Layered multicast and fair bandwidth allocation and packet prioritization

ABSTRACT

Embodiments include an overlay multicast network. The overlay multicast network may provide a set of features to ensure reliable and timely arrival of multicast data. The embodiments include a congestion control system that may prioritize designated layers of data within a data stream over other layers of the same data stream. Each data stream transmitted over the network may be given an equal share of the bandwidth. Addressing in routing tables maintained by routers may utilize summarized addressing based on the difference in location of the router and destination address. Summarization levels may be adjusted to minimize travel distances for packets in the network. Data from high priority data stream layers may also be retransmitted upon request from a destination machine to ensure reliable delivery of data.

CROSS-REFERENCE TO RELATED APPLICATIONS

This application is a divisional of pending U.S. patent application Ser.No. 13/619,969, filed Sep. 14, 2012, which is a continuation of U.S.patent application Ser. No. 12/486,656, filed Jun. 17, 2009, now U.S.Pat. No. 8,514,718, which is a divisional of U.S. patent applicationSer. No. 11/342,167, filed Jan. 26, 2006, now U.S. Pat. No. 7,733,868,which claims priority to U.S. Provisional Application No. 60/647,601,filed Jan. 26, 2005, all of which are incorporated by reference as ifset forth in full herein.

FIELD OF THE INVENTION

The invention relates to network management. Specifically, the inventionrelates to the management of data packets to support multicasting.

BACKGROUND

Despite the versatility in digital communication, interoperability andinternationally accepted communication protocols of the Internet, itsfundamental design has not changed much since its conception and doesnot excel in everything. Watching live TV for example is something whichis not typically done over the Internet, even though television has beenaround almost twice as long as the Internet Protocol and represents ahuge market. The reasons for this are based on the design of theInternet and Internet Protocol (IP).

The Internet is a packet-switching network where data is exchanged insmall units or packets that are independently transported over thenetwork and concatenated again at the receiver into its original form. Astrength of packet-switching is that it allows for very flexible use ofthe physical network wires. When two communicating parties have no datato exchange for a certain period of time, no packets are sent and thewires can carry packets from other parties. On the Internet, bandwidthis not reserved, but available to and shared by everyone. Theconsequence is that it cannot guarantee a minimum amount of end-to-endbandwidth, making live video streams often appear jerky because framesare skipped due to congestion that delays or prevents delivery.

Even though with help from specialized protocols such as Distance VectorMulticase Routing Protocol (DVMRP) or Protocol Independent Multicast(PIM), the Internet Protocol allows for data packets to be multicast toa large number of receivers simultaneously, using this feature tosuccessfully realize a live video broadcast is a challenge. A videostream is transmitted at a fixed high rate and not all parts of thenetwork are likely to have sufficient bandwidth available to forward thestream.

When a bandwidth bottleneck is reached, the router discards the packetsthat cannot immediately be forwarded. This causes two problems. The datastream that is eventually received by one or more receivers further downthe network is corrupt and the congestion also has a negative impact oncommunication sessions of other nodes that communicate through thebottleneck router. The only way to avoid this problem using the InternetProtocol and standard multicast is to find a transmission rate that issupported by all parts of the network. However, since the network isavailable to anyone, this rate will continuously change. A transmissionrate is selected and the packet loss is accepted. However, when packetsare dropped randomly by overloaded routers the data stream will sufferpacket loss. If additional packets are sent through the bottleneckrouter, there is a larger chance that the router will choose one ofthese packets when ready to send another packet, implicitly rewardingheavy streams during congestion. This encourages sending redundant datathereby exacerbating the problem.

A more fundamental problem of flow control using the Internet Protocolis that slowing down the data may not be an option for certain types oflive data streams. However, packet loss is unavoidable using theInternet Protocol and while data types such as audio and video data canusually withstand some packet loss without becoming too corrupted toplay, this does not apply to all types of live data. Real-time financialdata, for example, will become useless and even dangerous to use ifrandom packets of trades are lost.

BRIEF DESCRIPTION OF THE DRAWINGS

Embodiments of the invention are illustrated by way of example and notby way of limitation in the figures of the accompanying drawings inwhich like references indicate similar elements. It should be noted thatreferences to “an” or “one” embodiment in this discussion are notnecessarily to the same embodiment, and such references mean at leastone.

FIG. 1 is a diagram of one embodiment of an overlay multicast network.

FIG. 2 is a diagram of one embodiment of the basic components of arouter daemon in the overlay multicast system.

FIG. 3A is diagram of one example embodiment of a layered multicastsystem network divided into logical clusters.

FIG. 3B is a diagram of a routing table for the network of FIG. 3A.

FIG. 4 is a diagram of an example network with hierarchical structure.

FIG. 5A is a diagram demonstrating summarization of an example networkwith an S2 domain.

FIG. 5B is a diagram demonstrating the stretch factor for the examplenetwork of FIG. 5A and shows the inner network of the S2 domain.

FIG. 6 is a diagram of an example embodiment of a multicast distributionor “sink” tree.

FIG. 7 is a diagram of an overlay multicast network with a failed link.

FIG. 8A is a flowchart of one embodiment of a process or managingcongestion in the overlay multicast system.

FIG. 8B is a flowchart of one embodiment of a process for handling layerrepair.

FIG. 9A is a diagram of two router daemons connected by a link.

FIG. 9B is a diagram of two router daemons connected by a link where aretransmission is requested.

FIG. 10 is a diagram that shows the process of selective packet repairand ordering.

FIG. 11A is a diagram showing thinning of a stream of data.

FIG. 11B is a diagram of a multi-level token ring structure.

FIG. 11C is a diagram of a multi-level token ring structure showingthinning of data.

DETAILED DESCRIPTION

To provide multicast services a network needs to support one-to-manycommunication that can send data packets from a data source to more thanone receiver, ideally without putting extra stress on the network orsource when the number of receivers increases. Multicast routing can beoffered by different methods. One method is to let receivers tell thenetwork, but not necessarily the source, which data streams they want toreceive and let the network compute data distribution paths that deliverjust the right packets to each receiver. Multicasting can also be doneby letting the source encode the list of receivers in each data packet,thereby freeing the network from the potentially computationallyintensive task of maintaining multicast distribution paths. However,this method does not scale to handle a large number of receivers. Athird method relies on logic at the receivers by letting the networkapply a broadcast mechanism whereby each packet is delivered to everyconnected node and letting the receivers filter out only those packetsthat are interesting. This method may also generate a heavy load on alarger network but it is simple.

In one embodiment, a multicast network is constructed as an overlaynetwork. In one embodiment, the overlay network includes a number ofsoftware implemented routers connected by normal IP or TCP connections.A mesh is created in which every software router is connected to one ormore other software routers by means of virtual connections that appearto be direct connections to the other software routers, but are likelyimplemented by a number of intervening traditional TCP/IP routerssituated between the software routers. Two routers that are connectedthis way, are adjacent in the perspective of the overlay, but in realityare many physical hops away from one another. Also, a software routerthat has three software router neighbors has three independent virtuallinks. However, it is possible that this software router only has asingle physical network connection that is shared by the three virtuallinks.

To the underlying network, the overlay network is nothing more than acollection of applications that send data between static pairs. Beyondeach router pair (identifiable by their TCP connection) there is norelation between the individual daemons. As such, the overlay networkcan easily work through firewalls, NAT (IP masquerading), proxies andVPN's. Firewalls cannot control which software router POPs can talk toeach other. The system shares some principals with a HTTP proxy servertunneling traffic to and from web browsers. In an example HTTP proxyserver system, an intranet has two web browser machines. They bothbrowse the Internet using the proxy server also on the local intranet.Although the individual web browser machines can surf the netindividually, they never make a direct connection with a remotewebserver, but only with the local proxy server that acts like a relaypoint. As such, the firewall that sits between the proxy server and theremote webserver can only choose to either allow the proxy server totalk to the webserver or deny it. It is generally unable to enforceunique policies for individual browser machines. It generally cannottell on behalf of which client the proxy server is fetching a webpage.Returning to the overlay network, because each software router is arelay point that tunnels data traffic for different senders andreceivers similar to a HTTP proxy server, firewalls have no fine-grainedcontrol over communication over the overlay network. As soon as afirewall allows for only a single TCP connection between an internal andan external POP, all software routers connected to the internal one cantalk to all POPs connected to the external one, and vice versa, withoutrestriction.

In one embodiment, the overlay multicast routing system also managesflow control and timely delivery. Non interactive live data streams donot actively anticipate network congestion. To manage congestion thenetwork manages the available bandwidth to allow for fair or equaldivision among the streams. Without management, high volume streams areassigned a larger capacity percentage on overloaded links resulting inlittle benefit in keeping the bandwidth requirements of a stream low, asthe packet loss percentage is determined by how much the network isoverloaded by all streams combined and not by the requirements of theindividual streams. An alternative to letting the network handle theflow control and congestion is to put the responsibility at the sourceand receivers. However, letting data streams anticipate networkconditions requires a form of feedback information from the network orthe receivers. In this case it is beneficial that the amount of feedbackdoes not grow linearly with the size of the audience, as that wouldreduce the scalability of the multicast. Even when a scalable form offeedback information can be realized and the data stream adapts itstransmission rate according to the network conditions, the problemremains that live streams lose their value when they are slowed down anddelivered late.

In one embodiment, an overlay multicast system also implements ormanages delivery of packets including an option for guaranteed delivery.It would be ideal if every receiver would receive the data streamwithout loss or corruption. However, when the content is ‘live’ andcannot be slowed down, but the network has insufficient capacity, packetloss is difficult to avoid. In fact, even when there is sufficientbandwidth at all times, store-and-forward packet-switched networks arenot able to guarantee the delivery of all packets. For example, when arouter crashes, all packets in its buffers may be irrecoverably lost. Ifa network uses dynamic routing, packets are also dropped when routingpolicies change or packets are trapped in an occasional, temporaryrouting loop. In cases where packets are ‘accidentally’ lost, anend-to-end mechanism of retransmissions can be applied that cancompensate for the loss. However, since this requires a form of feedbackinformation, it is beneficial for reasons of scalability that theoverhead involved with retransmissions is not linearly related to thesize of the audience.

End-to-end retransmission feedback may be avoided in at least two ways.First, it is possible to let the network components keep a copy of themost recently forwarded packets and let them participate inretransmissions by intercepting the retransmission requests andservicing them locally. This approach often utilizes greater storage andincreased processing power requirements at the network components.

The second alternative to end-to-end retransmission requests is that ofencoding redundant information in the data packets. If enough redundancyis encoded, a lost packet's content may be entirely recovered from theextra information in the other packets. The downside of this system isthat it comes with a constant level of bandwidth overhead that isrelated to the level of packet loss tolerance, regardless of whetherpackets are actually lost. Each of these approaches to packet loss failin the case of a live data stream that produces more bytes than thenetwork can forward. When local or end-to-end retransmission requestsare used, the problem may even be increased as the retransmissionrequests use extra bandwidth, causing more data packets to be lost.

Embodiments of the overlay multicast system are designed to offer anend-to-end solution to efficiently multicast live data including stockmarket data to any machine connected to the Internet. The system iscapable of giving certain hard guarantees over what is delivered toreceivers. If packets must be dropped due to congestion or otherirrecoverable problems, it is done in a fully deterministic way thatdoes not corrupt the data. Where a receiver of data such as a viewer ofa film may accept the random loss of one or two video frames, this typeof data loss may wreak havoc in financial data when the missed packetcontains an important trade. The system supports deterministicallydelivery designated parts of a data stream when the network lackssufficient capacity. The system utilizes a layered multicast with datastreams subdivided into individual streams or layers. This allowsreceivers to subscribe to only those layers that the network can handle,so that random packet loss can largely be avoided.

In another embodiment, an enhanced form of layered multicast is usedthat guarantees complete delivery for certain layers to avoid randomloss altogether, making it suitable for certain types of critical datasuch as market data. The system is characterized as controlling twoprimary activities. The first activity is running and managing a robustand scalable overlay network that uses its own routing algorithms andsupports multicast, while the second activity is managing flow controland congestion when live streams overload the network and ensuringlayered multicast can be offered with guarantees.

FIG. 1 is a diagram of one example embodiment of an overlay multicastsystem. In one embodiment, the system includes an overlay network 100that includes a number of software router daemons 109, 115, that areinterconnected by normal TCP connections 119 or similar reliablecommunication connections. The overlay multicast system forms anintricate web of routers and virtual links similar to the Internet, butoperating on the application layer. This system operating at theapplication layer ‘overlays’ the physical network and other layers ofthe Internet providing its own system of routing, controlling the lowerlevels of the network. Any number of routers and client applications maybe a part of the system. In the example, two of the routers 109, 115 arein communication with local client applications 101, 111.

In one embodiment, each of the routers have a unique, string-basedaddresses 117. In another embodiment, the routers may have other typesof unique addresses such as numerical addresses. Each of the routerdaemons 109, 115 executes a routing algorithm that computes the shortestpaths that allows each router to send data packets to any other routerin the overlay multicast system as well as any machine in communicationwith these routers.

In one embodiment, the system includes a runtime client library 103,113, application programming interface (API) or similar system forsending and receiving packets that can be utilized by user applicationsto communicate over the system. This library connects an application toa nearby system router daemon 109, 115 through a TCP connection 107,121, native inter-process communication or similar communication method.The native inter-process communication method may be used if both therouter and the client application run on the same physical machine. Whenconnected to a router daemon, the client application 101, 111 can sendand receive data packets from the network under the router's uniqueaddress. Router daemons have a unique address, while user applicationsconnected to a router daemon are identified through logical ports.

In one embodiment, the topology of the overlay network may be configuredat router daemon startup. The topology may be relatively static afterconfiguration. In one embodiment, the relationship of router daemons toapplications may be one to many, with a single router daemon servingmultiple applications. In another embodiment, the relationship may bemany to many.

In one embodiment, the data packets in the overlay network may beaddressed to a single communication endpoint or to abstract multicastaddresses. Single communication endpoints are network addresses used byuser applications. Access to them is exclusive such that only one userapplication can use them for sending and receiving packets at a time.

In one embodiment, a unique network address in the overlay network thatis used by an application is the combination of a logical node addressassigned to the overlay router and the port name chosen by theapplication. If an application that is connected to an overlay routerwith the logical node address “n1.msx.1” wants to use a port named“myport” for receiving unicast data packets, the fully qualified networkaddress becomes “n1.mxs.1:myport.” Other applications connected to theoverlay network that want to send packets to this application can usethis as a destination address. Ports may be referred to as tports.Overlay network addresses must be bound, prior to sending or receivingpackets. When a data packet is sent from an endpoint address it willcontain this information as its source, allowing other routers or thereceiving application to send a response.

In one embodiment, the overlay network may support packets being sent toabstract multipoint destinations, or multicast addresses. An overlaynetwork multicast address is a destination that each machine on thenetwork may subscribe to. The overlay network software routers ensurethat a copy of each data packet published to this multicast address isdelivered to every user application that subscribed to it. In oneembodiment, multipoint destinations are single sourced. Only oneapplication can publish data packets to the multicast address, makingthe overlay network suitable for one to many, but not many to manycommunications. Because a tport session of a multicast address can befreely chosen, each multicast address explicitly contains the locationof the source on the overlay network. This makes multicast communicationless flexible because publishing is not anonymous, but it greatlysimplifies subscription management. In another embodiment, multipleapplications may publish to a multicast address and many to manycommunication is supported.

In one embodiment, both unicast and multicast addresses aresyntactically equal. For example, the address “n1.mxs.1:timeservice”could be a unicast address used and bound by an application thatprovides the local date and time in response to any data packet itreceives. However, it could also be a multicast address that anyone cansubscribe to. For example, subscription to the address may provideperiodic date and time broadcasts from the user application that boundthis address as a multicast address. In one embodiment, each packet maycontain a flag or similar indicator that indicates whether itsdestination address should be interpreted as a unicast or multicastaddress.

In one embodiment, the overlay network provides a set of higher layeredprotocols as services over the packet oriented base layer. Theseservices may be referred to as protocol endpoints. Any number ofprotocol endpoints may be defined and supported. In one embodiment, fiveprotocol endpoints may be supported. A unreliable unicast protocol(UUP), a reliable unicast protocol (RUP), unreliable multicast protocol(UMP), ordered layered multicast protocol (OLMP) and reliable orderedlayered multicast protocol (ROLMP) may be supported. The UUP offers abest effort unicast datagram service to applications. The RUP offersreliable unicast communication between peers. The UMP offers best effortmulticast datagram services to applications on the overlay network. TheOLMP offers multicast communication with receiver driven rate control.Complete delivery is not guaranteed, but the packets that are receivedare guaranteed to be in their original order. The ROLMP offers reliablemulticast communication with receiver driven rate control. Streamlayering allows each subscriber to receive the data stream in thehighest possible quality, while the source never has to slow down.

FIG. 2 is a diagram of one embodiment of the basic components of arouter daemon in the overlay multicast system. Each overlay networkrouter 109 includes a packet switching core or kernel 201. In oneembodiment, packets that are received, either from a connection to aneighbor router or from a connected application, pass through the kernel201. The kernel 201 may not handle some specialized control packets. Thetask of the kernel 201 is to inspect the destination of each packet anduse a set of routing tables to determine how to forward the packet.

In one embodiment, the kernel 201 forwards packets to neighbors througha set of interface modules 203, 205. In one embodiment, the kernel 201may execute on its own thread and be event driven. The kernel 201remains idle until it is notified by an interface module 203, 205 of anincoming packet or by an application 227 that is sending a packet. Thekernel thread is woken up, reads the packet from the interface module orapplication and processes it. If the kernel decides that the packet mustbe sent out through an interface, it passes it to that interface andwaits for the next notification.

In one embodiment, each router daemon in the network is connected to oneor more neighbors. This is done by establishing connections between therouters. In one embodiment, the connections may be long lived TCPconnections or similar communication connections. In one embodiment, arouter daemon 109 runs one interface module 203, 205 instance for eachconfigured neighbor or configured communication channel. In anotherembodiment, the router 109 may run multiple interface modules for aconfigured neighbor or communication channel or a single interfacemodule for multiple configured neighbors or communication channels. Theresponsibility of an interface module 203, 205 is to establish aconnection, e.g., a TCP connection, and to pass packets from the kernel201 to the connection and vice versa. Packets of the first kind arereferred to as outbound packets, while the latter are inbound packets.

In one embodiment, the kernel 201 maintains a unicast routing table thatis used for packet switching. To make it possible for the network tofind shortest paths as well as adjusting these paths when the underlyingphysical network's characteristics change, each interface module maymeasure the quality of its virtual connection. These measurements arepassed on to the kernel 201 when the link quality is found to havechanged. Inside the kernel 201, the measurements may be fed to a routingalgorithm to determine if the changed link quality alters any of theentries in the routing table. If the routing table is changed, it isadvertised to the neighbors by encoding the entire table or a portion ofthe table in a data packet and passing it to each interface. In oneembodiment, if this type of packet is received from a neighbor andpropagated to the kernel 201 through the receiving interface, the kernelinspects the message and analyzes it through the routing algorithm. Ifthe new information leads to changes in routing entries, the routersends its own related routing entries to all its neighbors.

In one embodiment, if a neighbor router crashes, the interface detectsthis through an error on the virtual link and passes an indicator to therouting algorithm. The routing algorithm then changes the costindication associated with the link to an infinite value or similarvalue indicating that the link should not be utilized. In oneembodiment, the kernel 201 does not distinguish between a crashedneighbor and an infinitely slow link. The kernel 201 only detects thatthe link is not to be utilized when reading the entry in the routingtable.

In one example, when a source application connected to router s wants topublish a live data stream to multicast group s:mygroup, whereapplications on node p and q want to receive it, the source may firstbind the group in publish-mode. Binding a multicast group inpublish-mode means that the group is owned by the binding process. Onlythe owner of the group is able to publish to it. The receiving or “sink”applications connected to routers p and q now bind group s:mygroup insubscribe-mode. Binding a multicast group in subscribe-mode results inthe router node being connected to a virtual multicast distributiontree. The subscribers receive a copy of any packet published by thesource.

In one embodiment, multicast groups do not necessarily need to be boundin publish-mode first. Any application may subscribe to any multicastgroup at any time. If there is no source, there will be no data toreceive. Binding a multicast group either in publish-mode orsubscribe-mode are distinct operations. When the source also wants toreceive a copy of its own stream, it binds its group in subscribe-modeand uses the resulting tsocket to read the data back. Data packets carryboth the address of their source application, as well as the networkaddress of their destination.

For unicast packets, the router uses its unicast routing table to findthe preferred next hop, while multicast packets are routed according toa subscription list in a multicast subscription table. In oneembodiment, a unicast data packet contains the node address of thesource router, the tport of the source application, the address of thedestination router and the tport of the destination application. Amulticast data packet contains the multicast group it was published to,represented by the node address of the source router and the groupidentifier that was bound by the source application. It does not containany other addressing information.

In one embodiment, the overlay multicast system determines unicastrouting tables and multicast subscription tables for each router. Thesystem utilizes any type of routing algorithms or protocols to determinerouting tables. Protocols that are utilized include distance-vectorprotocols and link-state protocols. Link-state protocols take therelatively simple approach of propagating the state of the entirenetwork as a list of all known links in the network with their cost.Distance-vector protocols are based on the Bellman-Ford protocol. Theywork by letting every router periodically advertise its own routingtable to its neighbors. In one embodiment, the Extended Bellman-Fordprotocol, hereafter referred to as ExBF, is used as the basis of theoverlay multicast network. For sake of convenience, the embodiments ofoverlay multicast system are described that utilize the ExBF, however,those of ordinary skill in the art would understand that other routingalgorithms may also be utilized.

In the ExBF protocol, a slight increase in the amount of information iskept for every destination. Instead of storing just the distance to eachdestination for every neighbor, ExBF also stores the address of thepre-final node of each path. Hence, instead of storing the collection ofdistances {Di(j)} where j represents a destination, D the distancebetween this node and j, while i ranges over the node's neighbors, eachrouter also stores {Pi(j)} where P is the pre-final, or so-called‘head-of-path’ node in the shortest path to destination j. Now becausethe router knows the pre-final node of every destination, it canbacktrack the full path to any destination by recursively looking at thehead-of-path of a destination and treating it as a new destination.

In one embodiment, the interface modules 203, 205 implement a simplealgorithm to establish the communication connection to the neighbormachine or device. The interface module 203, 205 attempts to connect tothe configured neighbor router by actively trying to connect to itsnetwork address, e.g., its TCP/IP network address. If the connection isaborted with a connection refused or other error, it is assumed that theneighbor is not yet running and the interface module 203, 205 starts tolisten on its configured IP port so that the neighbor can connect to itwhen it is started up. The interface module 203, 205 waits for anincoming connection request for a brief period of time. After this timeperiod expires, the interface module 203, 205 returns to activelyconnecting to its neighbor. To avoid a situation where both neighborscontinue to switch between the listen and active states at the sametime, the duration of the listening state is influenced by a randomfactor. An advantage of allowing each neighbor to switch between theactive connection and listening states when establishing a connection isthat it allows the connection of routers even if one of the routers ison a masqueraded IP network. A router on a masqueraded IP network isunable to accept incoming connections such as TCP connections.

In one embodiment, interface modules 203, 205 can be configured atrouter deployment through configuration files. This will make the routerdaemon automatically create the specified interfaces at startup. If theconfigured neighbor routers are online, all long-lived connections willautomatically be established. It is also possible to add new neighborconnections and interface module instances dynamically at runtime. Thisway the overlay network topology can be changed flexibly and new routerscan be added to the network. If an interface module 203, 205 manages toestablish a connection with a neighbor router, the interface modules ofthe routers exchange overlay network node address information to informeach other of their presence. When the interface module 203, 205receives the node address of its neighbor, it passes this information,together with an announcement that a connection has been made to thekernel 201. This information allows the kernel routing algorithms tobuild dynamic routing tables.

In one embodiment, a role of the interface modules 203, 205 is toestablish the connection with a configured neighbor router and to senddata packets from the router kernel 201 to the neighbor connection andvice versa. The interface module 203, 205 incorporates a framework thatallows custom software plug-ins to influence the stream of packets thatflows between the network and the kernel 201. This mechanism is referredto as the interceptor pipeline 251, 257. In one embodiment, eachsoftware plug-in component that needs to control the packet flow is aclass that implements a simple programming interface. In anotherembodiment, the software plug-ins may have any implementation structureincluding objected oriented structures. This interface allowsinterceptor instances to be chained together, forming a packetprocessing pipeline. The contract of an interceptor is that it receivesa packet, applies it operations and then passes the modified packet tothe next interceptor in the chain.

In one embodiment, an interceptor pipeline 251, 257 sits between therouter switching core and the network connection. When the router kernel201 delivers a packet to the interface module 203, 205 for transmissionto the neighbor router, the interface module runs the packet through theinterceptor pipeline, giving the interceptors the chance to modify thepacket. Each packet that comes out of the pipeline is transmitted to theneighbor router.

In one embodiment, each interface module 203, 205 has two interceptorpipelines 251, 257. The first interceptor pipeline 257 is used toprocess outbound packets. The second interceptor pipeline 251 is usedfor inbound packets. These pipelines are independent of one another, theordering of the interceptors and the number of processing steps can bedifferent for inbound and outbound packets. Each pipeline can beconfigured uniquely. Interceptor pipelines may have any number ofinterceptors, provided they do not add too much latency.

In one embodiment, an example of an interceptor is one that filterspackets from a specific overlay network router. When this interceptor isimplemented as a manageable component that can be configured dynamicallyat runtime, it can be used to implement basic firewall functionality onthe overlay network. If it receives a packet that matches its rejectionpattern, it discards the packet by not passing it to the nextinterceptor in the pipeline. Another type of interceptor that may beused is a traffic monitor that counts the size of each packet thatpasses by and uses this to log traffic activity and compute bandwidthstatistics. This plug-in mechanism allows an overlay network router tobe extended with additional functionality without modifications to theunderlying software.

In one embodiment, the interceptor pipelines 251, 257 act as a packetbuffer between the router kernel 201 and the network. The interfacemodules 203, 205 temporarily buffer inbound and outbound packets. Thisensures that the kernel 201 is not blocked when sending a packet to aninterface. In one embodiment, the interface modules have a separatethread or set of threads that continuously dequeue and serialize packetsfrom the buffer and write them to the communication connection orenqueue into buffer received packets.

In one embodiment, the interceptor pipeline provides a temporary packetbuffer and offers inter-thread communication between the kernel threadand the interface thread. Interceptors may be divided into twocategories: interceptors that block and interceptors that returnimmediately. The first category is referred to as blocking orsynchronous interceptors. In one embodiment, to avoid situations where arouter kernel 201 is blocked for an arbitrarily long time, aninterceptor pipeline may contain one non-blocking interceptor. Anon-blocking interceptor guarantees immediate return of control to acaller by storing the packet in an internal buffer. The packets in thebuffer may be discarded if the buffer exceeds a certain threshold size.

In one embodiment, a maximum size limit is placed on the packet buffersto prevent them from exhausting the router's memory. Storing packetsbefore processing them means that their delivery will be delayed. Thelarger the buffer gets, the longer the packets are delayed. Because ofthis, the interceptor drops packets when the buffer reaches its maximumsize. In one embodiment, the system uses reliable TCP or similarconnections for transmitting packets between routers in which packetsare only dropped inside the buffer interceptors of the interfacemodules.

In one embodiment, the interceptor plays a role in making packet lossdeterministic. Packets are not explicitly dropped in any other part ofthe layered multicast system network, except the buffer interceptor inthe interface pipelines. However, that may not guarantee that a packetthat successfully makes it through all buffer interceptors of thenetwork's routers is delivered at its destination. Aside fromcontrollable packet loss, the network may also occasionally lose packetsin a non-deterministic way, for example when a router crashes withpending packets in its buffers, or when a connection between adjacentrouters is unexpectedly closed during transmission of a packet.

In one example, an inbound interceptor pipeline 251 may be structuredsuch that traffic throughput monitor interceptors 253 are positionedafter a buffer interceptor 255 and a firewall interceptor 265 may bepositioned before for the buffer interceptor 255. In an example outboundinterceptor pipeline 257, a throughput limiting interceptor 259 may bepositioned before a buffer interceptor 261 and a traffic monitorinterceptor 267 may be positioned after the buffer interceptor 261.

In one embodiment, the overlay network is accessible to applicationsthrough the use of the client-side programming library 221. This libraryconnects to a router daemon 109 at application startup and communicateswith it using remote procedure protocol or similar communicationprotocols. The communication between the library 221 and the routerdaemon 109 is driven by the application 227. The application 227 invokesfunctions in the router 109 through the library 221. In one embodiment,when the kernel 201 receives packets addressed to a tport that is boundby the application 227, it stores them until the client actively picksthem up. The application 227 through the library 221 continuously pollsthe router for packets. To minimize the overhead of the pollingmechanism, the router poll function does not return until there is atleast one packet delivered by the kernel 201. If more than one packet iswaiting, the poll function returns all waiting packets at the time it isinvoked. In another embodiment, the router may send an indication suchas an invoking a marshaled stub from the client, event notification orsimilar indicator to the application 227 through the library 221 toindicate the reception of a data packet for the application 227.

In one embodiment, a user application 227 uses a client library 221through instances of the overlay multicast system communication sockets225. If a user wants to be able to receive packets, they reserve anetwork address. In one embodiment, a network address in the overlaymulticast system is represented as a combination of the address of therouter and a unique port identifier, reserved for the socket of therouter. For sake of convenience, a port in the layered multicast systemwill be referred to as a tport and a socket as a tsocket. When a userapplication creates a tsocket for receiving normal unicast packets, thetsocket automatically binds a local tport at the router daemon through aremote call to the client IO multiplexor 207 at the router. In oneembodiment tports are bound exclusively and other clients may not use itconcurrently.

In one example embodiment, communication sockets 225 communicate with alocal input output (TO) multiplexor 229 that coordinates the handling ofcommunication between the sockets and the router kernel through theformation of RPC calls, native inter-process communication or similarsystems. A local IO multiplexor 229 utilizes an RPC stub 223 or similarprogram to communicate with the router via an RPC skeleton 219 andclient adaptor 209. A client IO multiplexor 207 at the router managesthe relay of these socket requests to the kernel 201.

In one embodiment, the router daemon 109 uses a packet buffer totemporarily store packets for each connected client application untilthey are picked up. An interceptor pipeline in the client adaptor 209 orsimilar process may be utilized for this buffering function.

To overcome the problem of growing routing tables, computation time andexcessive advertisement overhead, large networks can be partitioned intosmaller sub sections, connected by gateways. While the gateway routersmaintain routing information necessary to route packets to nodes inother network sections, nodes inside a section or domain only maintaininformation for those nodes inside the same domain. By substitutinglogical ranges of hosts in the routing table by one single condensedentry, the size of the routing table is reduced. This is called addresssummarization. This mechanism introduces a form of hierarchy that allowsthe network as a whole to grow far beyond the practical limits ofstandard distance-vector or link-state algorithms. The farther adestination host is away, the more efficient it can be condensedtogether with other remote hosts. The more summarization is applied,though, the less efficient the paths become on average. The factor bywhich the actual data paths on summarized networks differ from theoptimal paths, is known as the stretch-factor: the maximum ratio betweenthe length of a route computed by the routing algorithm and that of ashortest path connecting the same pair of nodes.

In one embodiment, the overlay multicast system takes a relativelystraightforward approach to address summarization. An administratordecides at deploy time which nodes form clusters and which clusters formaggregations of clusters. This is done by encoding hierarchy in thelayered multicast system node addresses using a dotted or similarnotation. Node addresses may be ASCII strings. In one embodiment, thestrings are limited to at most 127 characters. In one embodiment, only[a-z] and [0-9] are available. In another embodiment any characters,numbers of similar symbols may be utilized. Addressing may becase-sensitive or case-insensitive.

FIG. 3A is diagram of one example embodiment of a layered multicastsystem network divided into logical clusters. The example illustrates anetwork of eight nodes, divided into three clusters. Assigning nodes toclusters may be based on geographical properties, administrativeboundaries, wide area links and similar considerations. For example,nodes inside a corporate network are all assigned the same logicaldomain, whereas a network that connects nodes from differentcorporations, would usually assign a separate domain to each corporatenetwork. Another criterion for assigning nodes is that nodes that oftendisconnect and reconnect again later, possibly because the overlaymulticast system routers run on personal computers, are placed in asubdomain, to avoid the routing updates triggered by their changingstate to propagate far into the network.

Given the logical domains that cluster groups of nearby nodes, each nodecan treat domains other than its own as a single entity and use awildcard address that matches every host inside that domain. Thisreduces the size of the routing table, as well as the amount of routinginformation that needs to be exchanged between the nodes when thetopology changes inside a domain. For example, when a new host is addedto domain S 1, there is no need to propagate that information to theother domains, as they already have a wildcard entry that will match thenew address.

FIG. 3B is a diagram of a routing table for the network of FIG. 3A. Theexample routing table shows the effect of address summarization in thisnetwork on the routing table of node S2.C. The cost value that is listedin the third column represents the cost of the shortest path to thenearest node inside that domain. In the fourth routing entry, the numberin the cost column is the cost to reach S1.R from S2.C. If it is assumedthat both interdomain links S2.0-S1.R and S2.A-S1.R have the same weightor cost, S2.0 will route all traffic for domain S1 through neighborS1.R.

In one embodiment, when nodes exchange distance vectors, either becauseof a link change, or as a normal, periodic exchange, the receiving nodefirst summarizes the destination address of each distance vectorrelative to its own address. Summarization is done by matching thedestination address with its own address field-by-field and when a fielddoes not match, the rest of the address is substituted by a wildcard.For example, when S2.0 receives a distance-vector from S1.R thatcontains a path to destination S1.P, it is immediately summarized toS1.* upon arrival at node S2.C. This is because the first field differsfrom the first field of the local address and as such the remainingfields are replaced by a wildcard. This wildcard value is then fed tothe distance vector algorithm that checks whether the new cost orpath-length is shorter than the cost of the entry that was already inthe routing table. In the present example there already is an S1.*wildcard entry in the routing table that was derived from neighbordestination S1.R. Since the path to S1.P runs through S1.R, the path toS1.R will always be shorter than the path to S1.P, so the entry in therouting table will not be changed and no further routing updates will bepropagated to S2.C's neighbors.

In general, when a local address is 1.2.3.4 and an incomingdistance-vector advertises destination 1.2.2.2, it will be summarized to1.2.2.*. Destination 2.1.6 becomes 2.*, destination 1.2.3.4.5.6 becomes1.2.3.4.5.*, destination 1.2.3.5 stays 1.2.3.5 and destination 1.2.3.4.5also stays 1.2.3.4.5. Note that 1.2.3 and 1.2.3.* are two differentaddresses. The first only matches the exact address 1.2.3 while thesecond is a wildcard that matches everything that starts with 1.2.3 andhas at least 4 address fields. This includes 1.2.3.4 and 1.2.3.4.5.6,but not 1.2.3. If this mechanism of address summarization is used in anExBF implementation that carries tuples containing destination, cost andhead-of-path attributes in its vectors, then the destination address issummarized, as well as the head-of-path address.

FIG. 4 is a diagram of an example network with a hierarchical structure.When the example network first starts to converge using the ExBFalgorithm, node B.X advertises the following routing information toneighbor A.S: DV_(B,X,A,S): {(B.X, *, 0), (B.Y, B.X, 1), (B.Z, B.X, 1)}.When A.S receives the vectors, it summarizes its entries. What remainsis the single vector DV_(B,X,A,S): {(B.*, *, 0)}. The routing table ofnode A.S now contains RT_(A.S): {(A.S, *, *, 0), (A.R, A.R, A.S, 1),(B.*, B.*, A.S, 1)}. All addresses are summarized before processing.This includes neighbor addresses. The consequence of this is that when anode has more than one connection with a foreign domain, both neighboraddresses will be summarized into the same wildcard. This leads toambiguities and nondeterministic routing when this wildcard is listed asthe preferred hop in a routing entry, as it cannot identify a singleoutgoing link. This problem is solved by assigning a local identifier toeach link and using these numbers in the preferred hop column, ratherthan the addresses of the links peers. The routing update A.S sends toA.S.Y contains DV_(A.S,A.S.Y): {(A.S, *, 0), (A.R, A.S, 1), (B.*, A.S,1)} and leads to A.S.Y's routing table RT_(A.S.Y): {(A.S.Y, *, *, 0),(A.S.X, A.S.X, A.S.Y, 1), (A.S.Z, A.S.Z, A.S.Y, 1), (A.S, A.S, A.S.Y,1), (A.R, A.S, A.S, 2), (B.*, A.S, A.S, 2)}. When A.S.Y has finishedupdating its routing table, it advertises DV_(A.S,A.S.Y): {(A.S.Y, *,0), (A.S.X, A.S.Y, 1), (A.S.Z, A.S.Y, 1), (B.*, *, *), (A.S, *, *),(A.R, *, *)} back to neighbor A. S where the asterisk indicatesunreachable in the last three records to avoid long-lived loops. Theseloops were detected through the normal back trace mechanism of ExBF.

In one embodiment, summarization means that only a single path for arange of remote nodes is maintained. The consequence of this is thatpackets will not always be routed according to the real shortest pathbetween source and destination. To illustrate this effect in themulticast network, consider the example network of FIG. 4. Node S1.Rreceives the distance-vectors of both S2.A and S2.C. And aftersummarization learns that both neighbors offer a route for wildcardS2.*. If it is assumed that both interdomain links (S1.R-S2.A andS1.R-S2.C) have equals costs, then S1.R will choose S2.A to be thepreferred hop for S2.* because of the fact that its address logicallycomes first. When S1.R needs to forward a data packet to S2.B, it usesthe S2.* wildcard entry and forwards the packet to neighbor S2.A.Unfortunately this is not the shortest path to S2.B, as S2. A first hasto route the packet through S2.C. Instead, S1.R should have sent thepacket directly through neighbor S2.C. This ratio between the length ofthe actual path and the length of the optimal path between the endpointsare called the stretch factor.

In one embodiment, path stretching on the overlay multicast systemoccurs when packets are forwarded between nodes that are in differentlogical domains. An entire subnet is treated as a single node withseveral outgoing links. Because a virtual node often contains a largenumber of nodes, connected by its own internal network structure, it issometimes better to choose a different interdomain link when sendingpackets to the domain.

Since the overlay network runs its own adaptive routing algorithms, thecontent streams through the network are constantly rerouted to avoid thenetwork's hot spots and congestion. This can be particularly useful onwide area networks that are used for very different types ofapplications and hot spots are dynamic (i.e., moving around). Anotheradvantage of having custom routing algorithms is the freedom ofsubstituting them with others in the future.

In one embodiment, the software routers implement load balancing insidethe network. Traditionally routing algorithms seek for “best paths”through the network and send data streams over these. However, it issometimes much more desirable not to just send streams over this optimalpath, but to also select several sub-optimal paths and divide the streamover all of them. This also avoids the optimal path from gettingcongested when the stream requires more bandwidth than this single pathcan provide.

FIG. 5A is a diagram demonstrating summarization of an example networkwith an S2 domain. When node S1.P needs to forward a packet fordestination S2.F, it will send the packet directly down its interdomainlink to the S2 domain. In this example both interdomain links areassumed to have equal costs. FIG. 5B is a diagram demonstrating thestretch factor for the example network of FIG. 5A and shows the innernetwork of the S2 domain. FIG. 5B clearly shows that the stretch-factoris quite high for packets from S1.P to S2.F, as the optimal path runsthrough S1.R instead. Although the overlay multicast system addresssummarization technique cannot guarantee a maximum upper bound on thestretch-factor, it can manipulate the stretch-factor by changing thesummarization opacity.

By default, an address is summarized after the first field that differsfrom the local address. However, if that is changed to the second field,the overlay multicast system can look inside other domains for onelevel. A node with an address 1.2.3.4 will then summarize 2.3.4.5 into2.3.*, rather than 2.* and 1.3.4.5 into 1.3.4.*, rather than 1.3.*.Doing this at least at the border nodes in the overlay multicast networkthat have the interdomain links reduces the stretch-factor under certaincircumstances. As the overlay multicast network was designed to beadministrated by independent parties, administrators are free toexperiment with different summarization opacity levels withoutjeopardizing the other subnets or domains.

One advantage of the summarization of addresses in routing tables isthat the number of links that are present between nodes from differentlogical domains is irrelevant with respect to the number of routingentries in the routing table of a distance-vector protocol or similarprotocols. Since these tables only contain destinations with a singleforwarding strategy, the number of interdomain links does affect thesize of the routing tables. If y is identified to be the number ofentries in a node's routing table, x to be the total number of nodes inthe entire network, n to be the number of entities (nodes or nesteddomains) inside a domain and m to be the depth of the hierarchy. Fromthis it follows that the total number of nodes in the network can becalculated with:∀x=n^(m) where nϵIN∧n>1 and mϵIN∧m>0

The relation between domain depth, number of nodes and domain densitycan be expressed by the following formula for a network with a uniformtopology:

$y = {{( \frac{\ln(x)}{\ln(n)} )*( {n - 1} )} + 1}$

Given this formula of the address summarization's effectiveness, it canbe shown that routing entries in any network node remains under 100 evenwhen the network as a whole grows to well over 10 million nodes. It canalso be shown that the domain size n has no spectacular effect on thescalability. While a small value n yields a relatively deeply nestednetwork hierarchy, which implies more routing entries, it also meansthat each hierarchy level only contains a small number of entities. Alarge value n in a network of the same size yields a relatively flathierarchy with few levels, each level contains a large number ofentities. As a routing table can only contain a natural number ofentries, its number remains constant as nodes are added to domains thatare already captured by the existing wildcard entries.

An advantage of using a string based node addresses scheme is its largeaddress space and flexible naming. New subdomains may be added at anytime as there is no need for predefined or fixed hierarchical levels.All nodes inside a flat domain can be reconfigured into nestedsubdomains without the need to reconfigure routers or compute netmasks.Also, as address fields are not bound to a limited number of characters,they contain location or company names to make administration of thenetwork easier and more transparent. An example might be the addressn1.amsterdam.level3.marketxs.node4, or simply n1.ams.13.mxs.n3 to keepthe address relatively short. The drawback of these kinds ofunrestricted addresses is the amount of bytes they require. Becauseevery data packet always contains a source and a destination address,the amount of overhead per packet is excessive at times.

In one embodiment, the string based node address scheme is implementedby letting the first bytes of each address header specify the length ofthe address in bytes, followed by the address itself as a normal 8 bitASCII string with a maximum of, for example, 256 bytes. In onealternative, null-terminated strings are used.

In another embodiment, a sufficiently large address space is used anddivided into logical subsets, similar to IPv4 and IPv6 addresses. In oneembodiment, a 64 bit address space is used and divided into 8 bit fields(identical to IPv6) used to identify logical subnets. In thisembodiment, addresses might look like 120-23-61-201-43-146-128-132. Tomake distinct addresses, IPv4 uses the dot “.” for separation, IPv6 usesthe colon “:” and the present system uses a hyphen “-”. When eachaddress field represents exactly one logical subnet level, this schemeprovides a simple, but at the same time somewhat limited, way ofaddressing. It means that the complete network can address as manyindividual hosts as IPv6, while grouping them in at most 8 levels of subdomains, where each sub domain can contain a maximum of 256 hosts or subdomains. The sub domains can be utilized to indicate a network hierarchyin the topology, including a geographical or similarly based hierarchy.

In one embodiment, to reduce the amount of bandwidth wasted on thisoverhead, a mechanism is used for substituting the address strings forsmall 16 bit values or similar small sized value when transmitting apacket to a neighbor. Also, a translation table is kept at both sides ofthe connection that is used to put the original address string back inthe packet before passing it to the router's packet switching kernel.This optimization works on a per-link basis and may be totallyindependent from the communication with other neighbors. When it is usedon more than one interface, each interface maintains its own translationtable.

If a packet is transmitted to a neighbor, all address strings arestripped from the packet header. For each address string a 16 bit orsimilar value is generated and stored in the translation table. These 16bits or similar values replace the original addresses and before thepacket is transmitted, its protocol version field in the first header ofits wire-level representation is changed to a higher version. Thisversion number is reserved for optimized packets that would beunreadable for routers without support for address substitution. Toensure proper translation by the peer, the interface first sends aspecial packet that contains the new address substitutions. Such apacket is tagged with the higher protocol version number and containsone or more tuples of address strings and their substitute values. Inone embodiment, such a packet takes the form:

Type 0x1 is used for issuing new substitutions and type 0x2 is used toinvalidate an earlier or unknown substitution. The system attempts toensure that this (or similar) packet is received by the peer prior tothe data packet itself, which is something that cannot be guaranteed byevery transport layer. TCP or similar guaranteed communication methodmay be used for communication between neighbor nodes in the overlaymulticast network.

In one embodiment, configuring an interface for address substitution canbe done manually, or automatically. In the latter case the interfaceuses a special handshake packet that is part of the higher protocolversion. In one embodiment, the handshake packet takes the form:

The version field has the higher protocol version. The type for thehandshake packet is 0x0 and the length is 4. A peer transmits such a (orsimilar) packet when the connection is first established. When thepacket is echoed back by the peer, the interface knows that the neighboralso supports substitution and goes into substitution mode. After aconnection with a neighbor has been lost and is re-established, thehandshake is performed again as the peer could have been replaced bysoftware that does not support substitution.

In one embodiment, entries in the translation table all have a time-outor similar tracking mechanism to ensure timelines and accuracy ofentries. After this period the entry is removed from the table and anypacket still using the substituted value will not be processed and causethe receiving node to respond with a special packet containing thesubstituted values that weren't recognized, allowing the peer tosynchronize translation tables by sending a packet containing thesubstitution tuples. Since the invalidation of timed-out entries isautomatic, both peers use the same timeout or similar values in trackingmechanisms. In one embodiment, when the translation table is full,entries are removed according to a last recently used scheme. Theaddress substitution is implemented in a substitution interceptor thatis the lowest interceptor or a part of the lowest interceptor or otherinterceptor reconstructing packets immediately on arrival in aninterface module of a router. Substitution introduces additionalprocessing overhead and enabling it is optional for each node.

In one embodiment, the overlay multicast system may be used solely forrouting. The overlay multicast system offers multicast support overexisting IP networks. The overlay multicast system goes beyond offeringplain multicast functionality and additionally focuses on supportinglive multicast streams with limited jitter and delay through thecombination of overlay multicast, multicast congestion control andlayered multicast.

In the overlay multicast system the ExBF protocol or similar routingalgorithm implicitly contains information on which links in the networkare child links. For example, this is a by-product of the long-livedloop avoidance mechanism in ExBF. In ExBF, a node will inform itsneighbors about the fastest route it offers to a destination S, exceptthe neighbor that is in the path towards S. Instead, an infinitedistance is advertised to this neighbor. This information is used by arouter to conclude that each neighbor advertising that node S isunreachable, is actually using it as their parent node towards S andhence will immediately mark those links as child links in the shortestpath tree rooted at S.

Neighbors that do advertise reachable routes to destination S are notflagged as child links and therefore not used when forwarding multipointpackets from S. In another embodiment, this implicit information isavailable through any distance-vector protocol that uses poisonedreverse or similar techniques. Both methods are used to quickly delivera copy of any multipoint packet to all nodes in the network. Adisadvantage of both mechanisms however is that a copy of every packetis delivered to every node, regardless of whether or not that node isactually interested in multipoint packets from that source.

In one embodiment, the overlay multicast system is equipped with asparse forwarding mechanism. The overlay multicast system network isused for a single source multicast for high volume market data, videostreams or similar data to be sent to large numbers receivers. In thesetypes of uses it is not necessary that receivers themselves or othernodes be capable of publishing data packets to the multicast group.Instead, for security reasons it is better to know that only the realsource can publish to the group. Given these requirements, together withthe fact that multi-source protocols are more complex than theirsingle-source alternatives, in one embodiment, the overlay multicastsystem preferably uses a custom single-source protocol for multicastdistribution on the network. In another embodiment, the overlaymulticast system supports multiple sources. For sake of convenience anembodiment of a single source multicast network is discussed. One ofordinary skill in the art would understand that the principles ofoperation can also be applied to a multi-source application as well.

FIG. 6 is a diagram of an example embodiment of a multicast distributionor “sink” tree. In the example embodiment, the multicast system startswith a source q that is active, but in this example, a sparse-modeprotocol, with no receiver yet subscribed, no distribution tree exists.When the first receiver application p₀ subscribes, its router initiatesthe reverse computation of one branch of the sink tree rooted at sourceq. It does so by first marking that p₀ has a local subscriber for q:G bysetting the bit in LSp₀ [q,S] (LSp[ ] is a local array at node p thatkeeps track of all subscriptions of the local user applicationsconnected to router p) to true and then sending a join (e.g., a <join(q:G)>) packet to the preferred hop towards source q. This neighbor p₁receives the <join (q:G)> packet and marks the link with p₀ as a childlink for multicast group q:G. It then forwards the packet to the nexthop p₂ in the shortest path towards q. Eventually node p_(n) sends the<join (q:G)> packet to q. On receipt, q marks the link on which thepacket was received as a child link for its local multicast group G andstarts forwarding all packets addressed to G over the link to nodep_(n). The multipoint packets addressed to G are inserted into theoverlay multicast network at node q by the user application thatpreviously bound the local multicast group q:G in publish mode. Nodesp_(n), p_(n−1), . . . , p₁ all forward the packets to their neighborsfrom which a <join (q:G)> packet was received earlier. Node p₀ has nochild links for q:G, but does have LSp₀[q:G] set, so it delivers thepackets to the local application. When another node p₁ becomesinterested in q:G while LSp_(i)[q:G] is not already set, it sends a<join (q:G)> packet to its preferred hop towards q. Let p_(i+n) be inthe path from p_(i) to q that receives the <join (q:G)> packet andsuppose that p_(i+n) is also in the path between p₀ and q. In this case,p_(i+n) is already in the established part of the sink tree of q, so itdoes not need to forward the <join (q:G)> packet further to q. Instead,it only marks the link on which it received the packet as a child linkfor q:G. In general, a node u only forwards a <join (S:G)> packettowards S if it is not already subscribed to S:G (hence, LS_(u)[S:G] isnot set and the collection of child links for S:G is empty).

The example sparse-mode distribution tree generated when processessubscribe to a S:Q group originally equals a part or all of the optimalsink tree rooted at S. However, when the underlying unicast routingalgorithm detects changes in the network performance and updates some ofthe routing table entries, the optimal sink tree changes accordingly andmay no longer be matched by the multicast distribution tree, renderingthe tree as well as the multicast performance suboptimal. An extremeexample of this is when a link that is part of the multicastdistribution fails entirely.

FIG. 7 is a diagram of an example overlay multicast network with afailed link. In the example a node u currently in the distribution treefor q:G only has a single route to q via w. The link between u and wfails. The node u concludes that it can no longer be part of thedistribution tree (as u and q are no longer connected) and invalidatesall child link subscription information. In case LS_(u)[q:G] is set(i.e., the subscription table indicates that u is subscribed to group Gfrom node q), u will rejoin the distribution tree as soon as a new pathto q is found. In another embodiment, the algorithm sends a notificationto the locally subscribed application to indicate that u is no longerconnected to q.

In the example, it is assumed that all links have weight 1, while thelink between p and w has weight 100. The result is that the linkconnecting p and w is not in the optimal sink tree rooted at q. Sinceonly t and v have a local subscriber (LS_(t)[q:G] and LS_(v)[q:G] areset), the sparse distribution tree in this example equals the dottedarrows of FIG. 7.

In the example, nodes w, u, r and s all have only a single route to q.Node p has a path through both s and w, while s is the preferred hoptowards q. When the link between u and w fails, u becomes disconnectedfrom q and invalidates the subscription information for link u-r. Node uinforms neighbor r of this fact as part of an immediate distance-vectorexchange. This routing update implicitly tells node r that its join forq:G through u is no longer valid and that it should look for analternative neighbor to rejoin.

In the example, an explicit leave packet to u is unnecessary. Since r isalso left with no alternative path to q and has no locally subscribedapplications, it invalidates its q:G subscription and is no longer partof the distribution tree. Neighbors t and s are implicitly informedabout this through the ExBF routing update or similar routing updatesent by r. Eventually w, u, r and s all leave the distribution tree,while nodes t and v schedule a rejoin when a new path to q is found.Again, they inform their local subscriber applications about this. Whenp receives the routing update from s, it loses its preferred hop for qand switches to neighbor w. The new cost to q is now 100 plus distance(w, q) and neighbor s is informed. Upon receipt, s discovers the newpath to q and informs its neighbors r and v. Node v then sends the <join(q:G)> packet to s. Nodes s, p and w then construct the first part ofthe new, optimal distribution tree rooted at q. When the new pathreaches r and t, t sends a join packet to r and r to s, reconnecting allreceivers to the data stream.

In one embodiment, recovering from link failure or similar communicationerrors is divided in three steps. In the first step the underlying ExBFor similar routing protocol starts a wave that propagates a link failurethrough the network. Depending on the actual topology, this takes up toN−1 packets, where N represents the number of nodes in the network. Theworst case time complexity for recovering from a link failure is 3N whenExBF or similar algorithm is used as the underlying unicast routingalgorithm. A best case time complexity would be 2N.

The example discusses a scenario where the system recovers from afailure of a link that is part of the multicast distribution tree. Moreoften however, links will not fail completely, but rather fluctuate inquality, causing the underlying routing algorithm to reassign preferredneighbors. In this example case it is not a requirement that thedistribution tree is changed to reflect these changes, as it is notpartitioned. However, since the tree becomes suboptimal, it is changed.

In one embodiment, the decision whether or not to recompute adistribution tree after a routing table update is made according to thequality of the used path. When a node p in the tree detects that theneighbor in the optimal path towards the multicast source has a costthat is only marginally lower than the neighbor that is currently p'sparent in the actual distribution tree, the overhead of recomputationand the risk of packet loss outweighs the increased performance of thenew tree. Also, since the quality of each link is continuouslyre-evaluated, the updated sink tree may only be temporal. Twostraightforward solutions to this can be used. Either the algorithm setsa threshold on recomputation so that subscriptions are only moved to theneighbor with the currently optimal path towards the source if thedecrease in cost is at least a ratio x, where x>1. A larger value of xthen postpone tree adjustments until a substantial improvement can begained, while a smaller x makes the tree actively follow the changingcharacteristics of the underlying network.

Another solution is to postpone recomputation of the tree for a periodinversely proportional to the cost decrease ration x. The latter has theadvantage that the optimal distribution tree is always guaranteed to bereached in finite time since the last routing table change.

In one embodiment, to run the tree building protocol, two new messagesare introduced: the join message and the leave message. Both messagesuse a GroupData control message of the following format:

This message is used to send a list of subscriptions to a neighbor nodeor to cancel an aggregated list of subscriptions with a neighbor. Theaction field (9th byte) is used to distinguish between join and leave.The 10th and 11th byte indicate the number of (S, G) groups in themessage. The action field is either “join” (0) or “leave” (1). In analternative embodiment, this is extended to “stale”, “dead”, etc., toindicate the state of the publishing application.

Many multicast applications require some level of reliablecommunication. Examples of this are uploading files to several receiverssimultaneously or replicating web server caches. Without guaranteeddelivery, a multicast transport service has limited applicability. Inone embodiment, the overlay multicast network utilizes standard deliverycontrol based on moderating the sending of data based on the lowestbandwidth available on the route to a destination. However, thisdelivery system is unsuitable for certain applications.

One example is financial data distribution, it needs a transport servicethat can multicast live data without packet loss or corruption andwithout substantial end-to-end delay when parts of the network aretemporarily or permanently congested. While live market data cannottolerate random packet loss, it can be thinned. Under certaincircumstances it is acceptable to omit certain stock quote updates. Anexample is a desktop application that displays real-time updates forhundreds or thousands of financial instruments. If all quote updateswere delivered to this application, this would require a substantialamount of bandwidth and would cause the value of the more volatileinstruments to change faster than a human can read. The data may bethinned through a process that involves inspection of the individualstock quotes and encoding them in individual data packets, tagging eachwith a priority value. For this application, as long as all packetslabeled with the highest priority number are received, the partialstream can be considered intact. Additionally, when all updates of thesecond highest priority are also received, the quality of the partialstream is increased usually in the form of less latency.

In one embodiment of the overlay multicast system, priority numbers areassociated with data packets. The priority numbers represent a logicallayer inside a data stream. Data provided by a multicast application maybe in the form of a stream of data. This data is subdivided intocategories or priorities based on the nature of the data. When packetsin a data stream are labeled with priorities in the range 0 to 3, thestream is said to have 4 layers. Also, the convention is to treat 0 asthe highest priority and 3 as the lowest. Any other system ofidentifying priority levels may be utilized including alpha numericindicators or similar identifiers. If a stream only contains a singlelayer, all packets are labeled with priority 0.

To software router daemons, a priority value of a packet is relative tothe stream and becomes relevant when a decision to discard data at acongested router is made. The priority value has no meaning other thanas a criterion for discarding packets on congested links. When anoutgoing link of a router has insufficient bandwidth to transmit allpending packets, it forwards only those packets with a designatedpriority or the highest priority. This technique is also applied to anydata type, including audio/video data and financial data. Using thissystem, the packet priority numbers cannot be misused by sources to givetheir data packets a greater chance of prioritized transmission bygiving them the highest priority to gain advantage over other sources.Packet priorities are only compared between packets that are part of thesame data stream.

In one embodiment, knowing that routers will use the packet prioritynumbers when making forwarding selections on congested parts of thenetwork, a source carefully divides its data packets over differentlayers or priorities, in a way that a subset of the layers still containa useful, uncorrupted representation of the data. This system isespecially useful when multicasting a live data stream with a high datarate to a large number of receivers, scattered over a heterogeneous widearea network. Receivers that are on congested parts of the network willthen receive the highest priority parts of the data only. Thiseliminates the need to lower the publishing rate to match the slowestreceiver, while still being able to offer live, uncorrupted, thinneddata streams to clients suffering from insufficient bandwidth. Exampleapplications of the system include audio and video codecs that dividelive multimedia content over layers to enhance the user experience overwide area networks.

In one embodiment, incoming messages are sorted by unicast or multicastsender address. As described above, a sender address is the combinationof source router address and application session. An example of a sourceaddress is n1.mxs.office.erik:video.an2 which could be used by a userapplication broadcasting a video channel. In this case, only the sourceaddress is relevant, not the (uni- or multicast) destination address.Each incoming message is added to the queue that holds messages sent bythat particular sender. If this is the first message from a sender, anew queue is automatically created to store it.

FIG. 8A is a flowchart of one embodiment of a process for managingcongestion in the overlay multicast system. The congestion can bemanaged at each individual router. In one embodiment, the congestion ismanaged at the interface module level in each router. Each interfacemodule has an inbound pipeline and outbound pipeline discussed above forprocessing inbound and outbound data. Each pipeline buffers data that isawaiting further processing. However, if either pipeline is unable tokeep up with the pace of incoming data that needs to be processed somedata must dropped.

In one embodiment, data is received as a set of data streams at eachrouter (block 851). The data is then buffered in the inbound pipelinebuffer (block 853). The same process applies to outbound data that isreceived from the kernel by the outbound pipeline. This data is storedin the outbound buffer. After the data has been stored a check is madeof the inbound or outbound buffer to determine if it is full (block855). In one embodiment, data is stored in the data structures in thebuffers that organize the data packets into a set of queues. Each sourceaddress, data stream or layer in a data stream has a separate queue. Aqueue is sorted with highest priority and oldest data packets at thefront of the queue. If the buffer is full then a decision is made todrop a designated amount of data in the form of packets from the bufferto make room for incoming data packets (block 857).

In one embodiment, a queue is chosen randomly or in a round robin tohave data dropped. In another embodiment, the queue with the most datais chosen to have data dropped. In a further embodiment, a weightingfactor is calculated to determine which queue is selected to have datadropped. The weighting factor is based on the amount of data in a queue,size of packets in a queue and similar factors. The weighting factorcounteracts unfair distribution that is caused by selecting a queue by around robin, random or similar method of selection. Data streams withlarge packets are unfairly affected by other methods because adisproportionate amount of data is dropped in comparison with otherqueues with smaller packets. Queue selection is also influenced by thesize and the amount of data that is intended to be dropped. A largequeue that can drop close to the amount of data desired is weighted forselection.

In one embodiment, data is dropped if a total amount of data stored inall queues exceeds a threshold value. This threshold value is set by anadministrator or is a set value. These systems enforce the fairallotment of bandwidth between data streams. Also, this systemimplements the prioritization of logical layers by dropping lowerpriority level layers when congestion occurs. If the buffer is not fullthen the pipelines may continue to store data in the buffers (block851).

In one embodiment, while the queues grow in size as packets arereceived, a background thread constantly dequeues packets from thequeues and transmits them over the network.

In one embodiment, the system that manages in- and output for the queuesis divided in two parts. The first part is run by the background threadthat constantly dequeues packets from the queues and transmits them overthe network, while the second part is in charge of queuing new packetsthat are to be transmitted. The latter also implements the logic thatdefines when and which packet should be discarded (due to a bufferoverflow).

The dequeuing part of the system, when selecting a packet fortransmission, only looks at the first packet (most urgent) of eachqueue. To make sure each stream gets an equal share of bandwidth, ittakes the individual packet sizes into account. When all packets (notethat only the first packet of each queue is observed) have equal size,the dequeuing thread simply selects a random queue with equalprobability and dequeues one of its packets. Since each queue has thesame probability of being selected, each queue will deliver an equalamount of packets per time unit. Hence, each source will transmit anequal amount of bytes per second.

Since individual packets can have any size between 1 and some determinedmaximum number of bytes, their size must be considered when the transmitthread selects queues. Queues with many large packets should generallyhave a smaller probability of getting selected than queues with lots ofsmall packets in order to keep the bandwidth division fair.

In one embodiment, only the first packet of each queue is considered torepresent its queue and give it a probability of setting selected thatis reversely proportional to its size in bytes. In an example wherethere are three queues, P, Q and R where the first packet in P is 100bytes, the first packet from sender Q is 300 bytes, while the packetfrom R is 500 bytes, to calculate the selection probabilities, firsttheir respective selection weights are defined. For example, the weightof the packet in P is computed by dividing the total size of P, Q and Rby P's size: (100+300+500)/100=9. Q's weight is 3 and R's weight is 9/5.The weights are converted into selection probabilities by dividing themby the sum of all weights. This gives P a probability of9/(9+3+1.8)=0.65, Q a probability of 0.22 and R of 0.13. When a queuebecomes empty after dequeuing, it is removed.

In one embodiment, to keep the delay introduced by buffering packetsunder control, the second part of the system enforces a maximum totalqueue size. The sum of all packets in all queues may never exceed thisthreshold. When a new packet comes in, it is always accepted at first.If necessary a new queue is created for it, or it is added to itsdesignated, existing queue. After adding a new packet, the total size ofall queues is checked. If it is larger than the configured maximum, thealgorithm runs a removal round in which it first selects a queue andthen tells that queue to remove one packet. If the queues are still toobig after removal of one packet, the process of selecting a queue andremoving a packet is repeated iteratively until the total queue size issmaller than or equal to the configured maximum.

In one embodiment, shrinking the queues is always done after a packetwas added, never preemptively. The reason for this is to let the newpacket immediately participate in the removal selection process, ratherthan discarding it or making room for it at the expense of the otherqueues. Contrary to de-queuing packets, where the algorithm tries toselect the “most urgent” packet, we now need to select the “leasturgent” one. This packet is determined by looking at the size of theindividual queues (the largest queue should generally be shrunk to keepresource division among the streams fair) and which queue can match therequired size most accurately. This comes from the fact that individualpackets can differ greatly in size, so removing the “least urgent”packet from queue P could result in freeing 8 kilobytes, while removingthe “least urgent” packet from queue Q yields 40 bytes of space forexample. Now if the queues in total exceed the maximum size by only acouple of bytes, it seems logical to remove the small packet from Q.This policy comes with a consequence, namely that it implicitly promotesthe use of larger packets. After all, streams with lots of very smallpackets are more likely to accurately match the amount of buffer spacethat must be freed than stream queues with few very large packets. Thisproperty may encourage developers to use larger packets, increasing theefficiency and throughput of the network. In short, when selecting aqueue for shrinking, the system favors queues that will expunge theleast amount of bytes in order to match the total buffer capacity andare large compared to the others.

In one embodiment, before the selection can be made, the absolute weightfactor for each queue is computed. The weight factor combines thequeue's size and ability to accurately match the overcapacity of thebuffer by removing packets. The latter is exposed in v. It representsthe sum of the sizes of the packets that a queue must remove in order toeliminate the buffer's overcapacity. For example, if a queue P has 10packets of 100 bytes each while the buffer currently has an overcapacityof 170 bytes (suppose the maximum is 10000 bytes, while all queuescombined add up to 10170 bytes), P would need to remove at least 2packets (2*100 bytes). In this case, v_(p) is 200. The formula tocompute the absolute weight factor of queue m is defined:

$w_{m} = \frac{( {\sum\limits_{i = 0}^{n}v_{i}} ) \cdot S_{m}}{V_{m}}$In this formula, n represents the total number of queues, v representsthe number of bytes that would be removed if the queue was selected ands represents the total size of a particular queue. The formula capturesthe direct relation to queue size and the inverse relation to the amountof bytes that would be removed by the queue if it was selected.

When the absolute weight factors of all queues have been derived, theyare converted to weighted selection probabilities. This is done bydividing the individual weight values by the sum of all weights. Theselection probability of queue m is expressed in P_(m).

$P_{m} = \frac{w_{m}}{\sum\limits_{i = 0}^{n}w_{i}}$After the selection probabilities have been computed, a roulette-wheelalgorithm is run to select a queue and let that queue remove its vbytes. If v is smaller than the current overcapacity of the buffer intotal (note that this implies that the selected queue became empty andwas thus removed automatically), another selection round is done untilthe buffers shrink sufficiently. Note that the queue selection processis implemented as an atomic operation. During selection rounds no newpackets may be added to or removed from the queues.

In one embodiment, the overlay multicast system is a layered multicastin combination with a scalable, selective packet retransmissionmechanism to offer a service that can meet the demands of real-timefinancial data and similar applications. Selective packetretransmissions are crucial so that the overlay multicast system isguaranteed that subscribers will receive one or more layers that arecompletely intact. When an occasional packet in a high priority layer ismissed, it is repaired. However, the receiver may decide not to attemptto recover missing packets from low priority layers that are missing dueto network congestion. Trying to repair these layers would requireadditional bandwidth, resulting in additional packet loss. Yet, as moreend-to-end bandwidth becomes available, the receiver detects this andrepairs additional layers, so more intact layers are delivered to theapplication.

FIG. 8B is a flowchart of one embodiment of a process for handling layerrepair. In one embodiment, the overlay multicast system supports areceiver-side multicast socket that receives the raw packets from themulticast stream and is in charge of repairing damaged stream layersbefore forwarding them to the user application (block 801). This methodincludes a retransmission request packet. When the receiver detects amissing packet in a layer that should be repaired (block 803), it sendsa retransmission request for this packet towards the source (block 805).

In one embodiment, detecting packet loss is done in the conventional wayby sequencing each packet with an incrementing number. In anotherembodiment, each layer has its own sequence number, so all packets inone layer are sequenced independent of the other layers. This way,packet loss is detected in each layer, regardless of the conditions andpacket loss in other layers.

To avoid a cascade of retransmission requests when a packet is droppedclose to the source, a tree-based negative acknowledgment is used. Thistechnique provides localized repair without communication with thesource, because an intermediate router that has a copy of the requesteddata packet responds by sending it again and suppressing theretransmission request. When the next router closer in the distributiontree receives the retransmission request it checks its buffer todetermine if the requested packet is still stored therein (block 807).If the requested packet is there, it is retransmitted (block 809). Ifthe packet is not in the buffer then the retransmission request isforwarded to the next router in the distribution tree toward the source(block 811).

In one embodiment, when layers are dropped due to congestion, theprogramming API that is exposed to user applications notifies the userby means of an exception or special return value when it is readingpackets. These notifications may not be fatal and only serve to informthe user that the quality of the stream has changed.

In one embodiment, localized repair works by having each router daemonstore a window of transmitted packets. Packets are either stored for afixed period of time, after which they are discarded, or a fixed amountof buffer space is reserved to store the most recently forwardedpackets. Since the overlay multicast network by default use reliableconnections between router daemons, packets never get lost while theyare in transit between nodes. The only place where packets are purposelydiscarded is in the interface buffer interceptors. Thus, there is noneed to buffer packets that were actually transmitted over the reliableconnection, as those packets are guaranteed to have reached the neighbornode. In one embodiment, a router daemon copies packets to itsretransmission packet store before those packets reach a bufferinterceptor that discards packets.

The router's buffer interceptors copy packets to be stored forretransmission. In one embodiment, localized packet repair in theoverlay multicast system are implemented by adding an interceptor toboth interceptor pipelines of each interface and letting thoseinterceptors simply store a copy of each packet. Additionally, when aretransmission request passes through the router on its way to themulticast source, the interceptor inspects the request packet and if ithas the requested packet in its temporal packet store, it does notforward the retransmission request to the next interceptor. Therequested data packet is injected into the network again.

FIG. 9A is a diagram of two router daemons connected by a link. Thereare two interceptor pipelines 903, 905 inside each interface 907. Eachinterface contains both an outbound 903 and an inbound 905 interceptorpipeline. The latter processes packets received from the network, whilethe outbound pipeline processes packets just before they leave therouter daemon. In one embodiment, the pipelines contain two differentinterceptors; the CCI (Congestion Control Interceptor) 909 and the PRI(Packet Retransmission Interceptor) 911. In another embodiment, a routerdaemon contains additional interceptor instances per pipeline.

The CCI 909 implements bandwidth allocation rules and is responsible fordiscarding packets. During normal operation the source 913 publishes itsmulticast packet stream to its router daemon 901 (path 1) that sends thepackets to all interfaces that lead to interested receivers according tothe multicast routing table. In the present example, the packets arepassed to the first outbound interceptor (PRI) of the interface (path2). The PRI is responsible for storing a copy of each recoverable packetthat is transmitted. It stores the copies in the private temporarypacket store 915 of its interface (path 3). Then the interceptor passesthe packets to the next interceptor (CCI) (path 4) where it is buffereduntil the connection link with the adjacent router daemon has time totransfer it.

In one embodiment, if the packets are not dropped in the outbound CCI,they are received by the neighbor router daemon (paths 5 and 6) and sentthrough that router's inbound interceptor pipeline 921 (path 7). Theyfirst enter the PRI 923 (path 7), which stores copies in its inboundbuffer 924 (path 8) and then passes them on to the next interceptor(CCI) 925 (path 9) where they are temporarily parked until the router'skernel 927 thread picks them up (path 10) and delivers the packets tothe subscribed multicast tsocket which forwards them to their nextdestination 929 (path 11).

FIG. 9B is a diagram of two router daemons connected by a link where aretransmission is requested. If a packet is lost in the outbound packetinterceptor pipeline of the source router because the connection betweenthe routers was not fast enough to transmit all the packets, thereceiving tsocket notices the loss and send a retransmission request(path 1). The destination address of this unicast retransmission requestpacket is that of the multicast group. The network packet switchingkernels recognize the retransmission packets and use the multicast groupto route the packet towards the source of the group. The retransmissionrequest packet of the example passes the router kernel 927 and reachesthe outbound packet interceptor pipeline 931 and goes into the PRI 933(path 2). This interceptor scans specifically for retransmissionrequests and tries to answer them locally. To this end it inspects thepacket, looks up the multicast group, the stream layer identifier andthe layer's sequence number and then checks the inbound packet buffer935 to see if it contains the packet (path 3). If this is not the case,the packet is passed on to the rest of the interceptors (path 4) andeventually transmitted to the adjacent router (paths 5 and 6) where itis fed to the inbound interceptor pipeline 941 (path 7).

In one embodiment, the PRI 943 checks the interface's outbound packetbuffer 945 (path 8) and finds the multicast packet that was dropped bythe congested outbound CCI 909 earlier. It is then re-inserted into theoutbound packet stream via the outbound PRI 911 (path 9). Theretransmission request packet is then dropped and not forwarded anyfurther. Assuming the packet is not dropped again, it travels the normalway towards the receiver (paths 10, 11, 12, 13). In one embodiment, thereceiver's PRI sees the multicast packet for the first time and stores acopy in its outbound packet buffer (path 14). The packet then continueson to destination 929 (paths 15, 16, 17). This mechanism of unicastretransmission packets and interceptors that offer localized repair isflexible in the way that not all router daemons need to support localpacket repair. A router that does not support it simply forwards therequests and lets the upstream routers handle them. The network as awhole becomes more efficient when more routers support it.

In one embodiment, because storing a copy of every multicast data packetrequires storage capacity, the overlay multicast system links the packetbuffers of the packet retransmission interceptors to a central datastructure that will only store unique packets. This is done because inmulticast transmissions a router that is a branch point in the multicastdistribution tree otherwise ends up storing at least three copies ofevery data packet. This is because the packet is received on oneinterface and sent out over at least two other interfaces. If everyinterface would individually store its packets, the inbound packetbuffer of the receiving interface as well as the outbound packet buffersof the forwarding interfaces contain the same packets.

In one embodiment, the central packet store eliminates duplicate packetsby storing each packet only once and by keeping reference tables thatpoint to packets for every interface. When an interface buffer stores anew packet in the central store, while that same packet was alreadystored by another interface, the central packet store merely adds apointer to that packet instance to the interface's packet referencetable.

In one embodiment, each entry in the reference tables has a time-outattached to it. This is to ensure packets are only temporarily stored.In one embodiment, the same packet could pass different interfaces atdifferent times. The time-outs are not attached to the packets in thecentral store but rather to the packet references in the interfacereference tables. In one embodiment, packets are only expunged from thecentral store when all references to the packet have timed out. Anotheroptimization that can substantially reduce the amount of requiredstorage space is to only store packets that were explicitly discarded bythe Congestion Control Interceptors (CCIs) because of bandwidthconstraints. Storing other packets has little value, as those areguaranteed to have arrived at the next router. No retransmissions may berequested for them unless some non-deterministic packet loss occurs as aresult of a crashing node.

In one embodiment, if a packet was lost close to the multicast sourceand because of this and all receivers simultaneously send aretransmission request, the upstream routers still apply conventionalnegative acknowledgement “nack” suppression techniques to combine allconcurrent retransmission requests from its child links into a singlerequest that is forwarded upstream towards the source. Applying localpacket stores at every router reduces the latency of packet recovery.

In one embodiment, time to live values are adjustable per router, setsystem wide or similarly configured. A short time-out minimizes storagerequirements, but will also lead to more overhead towards the source asmore retransmission requests need to be propagated further upstream.Large timeouts will be more tolerant to late retransmissions, butrequire more storage. When a time-out of 60 seconds is used by therouters that means a lost packet can still be recovered after a minute.However, when the data stream contains real-time data such as stockquotes, such a delay in delivery cannot be tolerated. Problems get worseif the data packets also have to be delivered to the user applicationsin their original order. In that case the user does not receive any datafor up to a minute while the missed packet is recovered and all newerpackets are waiting for it. In one embodiment, a ten second timeoutsetting is utilized.

If all routers in the network use the same time-out value, there is nouse in propagating these requests as the upstream routers are likely tohave purged their copies as well. To avoid nack-implosions from slowreceivers right after a packet is purged, each router interfaceremembers which packets it has purged. When a retransmission request isreceived for such a packet, the interface may respond by sending aspecial packet to the receiver that indicates that the requested packethad timed-out and cannot be recovered anymore. The retransmissionrequest in this scenario is not propagated further upstream. When areceiver gets this timeout notification, it knows it will not have towait any longer for the missed packet. It will notify the userapplication of the fact that data was lost and continue to deliver thenext packets. Whether or not an application can tolerate packet losswill depend on the type of content. Audio and video will usually not beseverely impacted by the loss of an occasional packet, while real-timestock quotes become dangerous to use when it is not known what quote waslost.

In one embodiment, where layered multicast is combined with localizedpacket retransmission, a service is realized that applies packetrecovery only to those layers that can be delivered with the currentbandwidth capacity. When the network has insufficient resources todeliver all layers of a stream, it is beneficial if the receivers knowthis and will not attempt to repair all missed packets from all layers.Instead, the receivers may use a mechanism that provides them feedbackabout the current network capacity and use that to decide for whichlayers it will recover missed packets. The algorithm monitors the statusof the total stream and from this information derives how many layers ofthe stream can be reliably delivered to the user without stressing thenetwork. It then marks these layers as being intact, reliable orsimilarly labeled and sends retransmission requests when an occasionalpacket is lost from these layers.

In one embodiment, the algorithm only delivers packets from the reliablelayers to the application, but not before they have been put back intotheir original global order. In order for the algorithm to decide whichlayers are safely marked as intact, it requires some status informationabout the reception of the stream as a whole as well as the networkconditions. To provide this, every packet contains the number of themost recent packet from all other layers or similar sequence data. Asidefrom carrying its own sequence number, each packet contains the currentsequence number of all the other layers as well. By inspecting thesesequence numbers each time a packet arrives, the receiver determineswhether it has missed any packets from the layers it marked as reliable.If that is the case, these packets may be recovered. Each time a missedpacket is detected, a countdown timer is started for it. If the packetis recovered before the time-out expired, the timer is canceled andremoved.

If however the timer manages to expire, the layer is consideredimpossible to repair and may be, together with all higher reliablelayers (lower in priority), removed from the list of reliable layers.How long the time-out interval should be can be application or contentdependent. The interval determines how long delay is tolerated by theuser. Setting the interval to a low value means a temporary congestionmay corrupt a layer long enough to cause the receivers to drop it. Sincepackets occasionally get lost when unicast routing tables converge,changing the shape of multicast distribution trees, or when a routercrashes that had pending packets in its interface buffers, appropriatetime-out values are adjustable by an administrator according to networkconditions to be found experimentally.

In an overlay multicast network that uses packet prioritization, routersexplicitly introduce out of order delivery when a burst of packetsqueues in the outbound interceptor pipeline of a router interface with aslow connection. The time-out may be large enough to allow for this.Setting the time-out to a long period makes the stream much moreresilient to congestion, but also increases the time for the receiver todiscover that a layer must be dropped due to bandwidth constraints.Until congestion is finally detected, delivery of the previouslyreceived packets that causally depend on missed packets is postponed.Starting timers when packet loss is detected allows for congestiondetection. However, it cannot be used to guarantee low end-to-endlatency.

When no packets are lost and all are received according to their globalorder, the receiver cannot measure the total transmission delay. Assuch, a receiver cannot distinguish between an idle and a crashedsource. In one embodiment, aside from removing layers from thereliable-list, the algorithm is also able to detect when more bandwidthbecomes available and new higher layers can be added to this list, sothat a higher quality stream can be delivered to the user. This is doneby passively or actively monitoring the layers that are not currently inthe reliable-list and not under repair. Every time a packet is receivedfrom these layers, this fact is stored for a period that is equal to therepair time-out discussed earlier. When a packet from a layer isreceived with sequence numbers that shows that a packet has been lostfrom a higher layer that is not in the reliable list, no retransmissionrequest is sent, but the fact that this packet should have been receivedis recorded and stored for the same repair time-out period.

If during the time-out period the packet is received after all, possiblybecause it was delayed by a router, the stored record is marked asreceived. Because the state of each packet from every layer is recorded,the overlay multicast system builds a packet arrival history that isused to determine whether the reception quality was high enough to add ahigher layer to the reliable-list and repair any further packet loss. Inthe overlay multicast system the reception history or similar data maybe used to calculate a moving average that describes the amount ofpacket loss over the last x seconds, where x is equal to the repairtime-out. With a small moving average history, the receiver will quicklyrespond to increased network capacity, while a longer history will onlyadd layers when the network's capacity was sufficiently high for alonger period of time. The receiver may keep track of the receptionquality of each layer by inspecting the list of sequence number that isattached to each data packet.

This causality information adds overhead to each packet, linearlyrelated to the number of layers used in the stream. For example, whensequence numbers are 32 bits and the publisher uses all 256 availablelayers, each packet comes with a kilobyte of causality information,which contributes to at least 12.5% for an 8 Kb packet that is filled tothe brim. To ease layer administration in the receiver socket, anoverlay multicast stream only supports a static number of layers. When apublisher binds a multicast group address for reliable, layeredcommunication, it may explicitly specify the number of layers that willbe used.

In one embodiment, by default, packets are delivered to the userapplication in their original order. Not only are the packets inside theindividual layers restored to their natural order using their sequencenumbers, but the total ordering across the layers may also be restored.When the source sends three packets with different priority (each in adifferent layer), all three are received by the user application in theexact same order. A prioritized, layered packet is described asP_(1 (9, 4, 6)) where 1 represents the packet's priority or layer, 9 thesequence number of the last packet of layer 0 (the highest prioritylayer) that was sent at the time this packet was published, 4 representsthe sequence number of this packet and 6 represents the sequence numberof the last packet of layer 2. Also, the packet tells that the streamuses 3 layers in total: layer 0 up to and including layer 2. It isconcluded that this packet causally depends on packet 9 from layer 0 andpacket 6 from layer 2 and will only be delivered to the user applicationafter those packets have been delivered. In another embodiment, theoverlay multicast system does not reorder the packets, but may leavethis task to the application using the network.

FIG. 10 is a diagram that shows one embodiment of a process of selectivepacket repair and ordering. For purposes of explanation the notation {0(9, 2, 6); 1(9, 3, 6), . . . } to describe a sequence of packets whereP_(0 (9, 2, 6)) was sent prior to P_(1 (9, 3, 6)), is used for sake ofconvenience. The first segment 1001, shows the sequence of packetsoriginally published by the source. It reads from left to right, soP_(0 (9, 2, 6)) was published first, followed by P_(1 (9, 3, 6)), etc.In this example the source has published eight packets, divided overthree priority levels or layers (0, 1 and 2, where 0 is the lowest layerwith the highest priority). The second segment 1003 shows the packetstream as received by one of the subscribers. It shows two lost packetsand an incorrect ordering. Before the packets are delivered to the userapplication, they are stored in internal buffers during the repair andreordering process. This state is depicted in segment 1005. Here thepackets are in separate buffers, each representing a layer. The packetsare ordered inside the buffers. The illustration shows the missingpackets P_(1 (9, 4, 7)) and P_(2 (9, 4, 8)).

If the receiver currently only has layers 0 and 1 in its reliable-list,it will attempt to repair the hole in layer 1 by sending aretransmission request. Note that if the given sequence was just asnapshot of a running stream, the receiver would have detected themissing packet P_(1 (9, 4, 7)) when P_(0 (10, 4, 8)) was received,because this packet says it depends on packet #4 from layer 1. So evenbefore P_(1 (10, 5, 9)) from layer 1 was received in our example, thereceiver already detected loss in layer 1 and immediately scheduled arepair timeout for the missing packet and sent a retransmission request.In fact, if it is assumed that the receiver had received packet #2 fromlayer 1 prior to our snapshot of FIG. 10, then the conclusion would bethat packet #4 as well as packet #3 were lost.

In the example, shortly after P_(0 (10, 4, 8)) was received, packetP_(1 (9, 3, 6)) is received. The receiver places the delayed packet inthe appropriate receiver buffer and cancels the repair time-out itstarted earlier when it detected that the packet was missing. This isillustrated in segment 1007. Note that the hole in layer 2 may also bedetected when P_(0 (10, 4, 8)) is received, as that packet claims to besent after packet #8 of layer 2 was published, so either packet #8 fromlayer 2 got lost in the network, or was delayed. However, since layer 2is not in the reliable-list, a retransmission request is not sent.However, as for all packets, a timer is started for packet #8 of layer2. When the user application is ready to read packets, the algorithmreturns only packets from layers that are in the reliable-list. Eventhough some of the layer 2 packets were received, they are discarded andnot delivered. The resulting stream of packets that is delivered to theuser is equal to the stream originally published by the source, with alllayer 2 packets removed.

In this example, despite the fact that the network has insufficientcapacity, having dropped packets from every layer and delivered thepackets out of order, the user received a deterministic subset of thestream that is uncorrupted. Because the overlay multicast system sourcesare live data streams that cannot slow down or pause and becausereal-time data are not buffered by the source too long, receivers willonly attempt to recover lost packets for a limited period of time.Whether a packet is received, recovered or lost, its buffer slot will befreed after this timeout, resulting in a natural size limit of thereceive buffer. How large the buffer can get is related to the averagespeed of the stream and the length of the repair time-out.

In one embodiment, the algorithm implementation does not enforce a hardsize limit during packet repairs. A more troublesome situation occurswhen the user application does not read data from the socket fastenough. When this happens, the amount of pending data builds up in thesocket and both storage requirements and the transmission delayincreases. In one embodiment, this may be handled by removing a layerfrom the reliable-list, causing those packets to be discarded from thebuffers, while decreasing the amount of data that is delivered to theapplication. If a clean feedback algorithm that can keep the number oflayers in balance with the application's reading speed is not used, theoverlay multicast system throws a fatal exception to the user orprovides a similar indicator to the user application and closes thesocket when the receive buffer reaches a certain maximum size, or whenthe total time between arrival of packets and their actual deliveryreaches a threshold.

Although restoring the original global packet ordering before deliveringthe data to the user application is assumed to be appropriate for mosttypes of data, there is content for which each packet invalidates allprevious packets. This is the case for stock quotes and similar timesensitive data. For example, when a new stock quote update is receivedfor a financial ticker symbol, it renders the previous updates useless.For most applications that process real-time financial data, only themost recent information is interesting. When global ordering isrestored, the algorithm will postpone the delivery of the most recentdata until all prior packets have been received also. For a typicalapplication, such as a market data terminal, that merely displays quoteupdates to the screen, this postponing adds little value. When the burstof pending updates is finally delivered by the tsocket, the applicationupdates the symbol's last value on the screen, thereby leaving only thelast and most recent update visible and overwriting all pending updatesimmediately.

In one embodiment, because the reordering process adds additional delayto the data delivery, it can be switched off by applications that do notbenefit from it. Without reordering, global (causal relations betweenpackets from different layers) and local ordering (order of individualpackets inside a single layer) is ignored and packets are delivered tothe application immediately after they have been received. Disablingreordering has no impact on the reliability. Lost or delayed packets maystill be recovered for all layers the reliable-list, only the recoveredpackets may be delivered with additional delay. Whether or not thismakes them useless is up to the application to decide.

For example, with market data it is useful to know whether an update isolder or newer than the one previously received. This is because a stockquote may be overwritten by a newer one, but not the other way around.In this case, the source could add a logical time stamp or similarsequence indicator to each stock quote, so the receiver can decide howto handle the update. In one embodiment, aside from configuring alayered multicast tsocket to restore global ordering or no ordering atall, a receiver can configure a tsocket to only restore local ordering.Whether or not this is useful will depend on the type of application andthe content, but it offers the advantage that delivery of packets fromlower (high priority) layers are not delayed during recovery of packetsfrom higher, lower priority layers.

In one embodiment, an example market data publishing application tracksmarket financial updates in a linked list or similar data structure.Every time a new update is received, it is stored in the appropriateslot, overwriting the previous update. This way, the linked list alwayscontains the most recent trade for every symbol. A virtual tokentraverses the list at a fixed speed. If the stream is to be thinned toone update per second, the token will visit one slot every second. Whenthe token visits a slot, it removes the quote update and forwards it,leaving an empty slot. Empty slots will be skipped by the token withoutany delay. When one of the symbols has two updates per second in theincoming stream, the second incoming update after the last visit of thetoken will overwrite the update pending in the slot and the olderpending update is dropped. This allows the leaky bucket algorithm toalways provide the thinned clients with the most recent, live update ofevery symbol. While the virtual token visits the slots that have arecent quote update pending to be sent out, a thread receives theincoming market data stream and inserts new trades in the slots. Tolimit the outgoing transmission rate to one quote per second, the tokenthread sleeps for one second after it moved and sent an update from aslot. When the token is able to complete one full circle between twoincoming updates of every symbol, there will be no data loss. Instead oflimiting the outgoing bandwidth to a fixed maximum, the dequeuing threadcould also be configured to run as fast as possible, relying on externalfeedback such as a blocking write call from the network, to reducepublishing speed. This could be particularly useful when the algorithmis used on the server side of a point-to-point TCP connection to aclient application. In one embodiment, the server uses the algorithm tosend the data at original speed, until network saturation slows down thelinks. A potential problem with this technique is that the originalorder of the updates is lost. In fact, when every symbol updates morethan once while the token circles the list, the token will find everyslot filled with a quote update, causing the order of the output streamto match the fixed order of the slots in the linked list. For quoteupdates from different symbols, this is usually not a problem, as theyare not related to each other.

In another embodiment, an alternative algorithm may be used thatprovides another way of thinning a stream of data, such as live marketdata, with varying bandwidth to a fixed maximum rate, without the riskof providing stale data, for example stale market symbols. Thisalgorithm may be used to produce enhancement layers that form a layereddata stream, such as a layered market data stream. One example method ofdoing this is to have several instances of the algorithm running thateach receive a copy of the input stream, letting their token run atdifferent speeds and tagging their updates with a layer number. Thelayers may then be combined into a single stream. Unfortunately, thismay introduce redundant information in the stream because the higherlayers contain updates that are also available in the lower layers. Thistechnique of encoding an entire stream in different qualities andsending them to clients concurrently to meet their individual bandwidthcapacities may be referred to as simulcasting.

In one embodiment, the thinning algorithm may be modified to have eachlayer contain only information that is not already present in the lowerlayers, thereby eliminating redundancy among the layers. To achieve thisthe leaky bucket algorithm may be extended to a multi-level ring asillustrated in FIG. 11B, creating a virtual cylinder of verticallystacked, circular lists, where each ring represents a layer of theoutput stream. Each ring has its own token. In one embodiment, thetokens all circulate their ring at the same speed. In other embodiments,the tokens in each ring may be set to circulate at different speeds.This data structure also identifies columns. A column is the collectionof vertically stacked slots that can all store an update for the samesymbol.

FIG. 11A is a diagram of an example method of thinning a stream of livemarket data. The gray slots represent symbols for which a recent quoteupdate was received, while the transparent slots are empty. FIG. 11B isa diagram illustrating the token ring data structure. Multiple instancesof the thinning algorithm and data structure may be conceived as beingvertically stacked, each instance represents a stream layer and eachsymbol has a column that can store more than one update. FIG. 11C is adiagram of one embodiment of a representation of the stacked thinningalgorithm of FIG. 11A at work. In the example, the gray slots contain apending quote update, while the transparent slots are empty. Theillustration shows how the updates of a volatile symbol are spread outover several higher layers.

In one embodiment, the thinning algorithm has a single input thread thatreceives the original stream. Newly received updates may be inserted inthe lowest ring. If a new update arrives for a slot in the lowest ringthat already contains an update for that symbol, this older update isvertically shifted into the next ring, so that the lower slot can alwaysstore the latest update. Once an update is shifted upwards, it cannotreturn to a lower ring anymore, even if all other slots beneath it wereemptied.

In one example embodiment, symbols that have more than one trade duringthe time it takes the token to complete a full rotation of the ring willbe spread out to two or more higher layers, while symbols that have veryinfrequent trades will only occupy a single slot in the bottom ring. Anexample result of this process is depicted in FIG. 11C. FIG. 11C showsthe thinning algorithm using three layers to thin a market data streamthat contains updates for eight symbols. Four symbols are included indata structure and assigned to the four visible columns in the front.The figure also shows that all four symbols have pending updates in thebottom ring, while the ‘MSFT’ symbol is so volatile that it shiftedupdates into all three layers before the tokens could dequeue any of thepending updates.

In a case where a symbol is so actively traded that it manages to shiftan update out of the top ring, that update is irrecoverably lost. Ifrequired, loss of updates can be avoided by letting the token in the topring run without dequeuing delay, letting it empty every slotimmediately after an update was queued. The result is that thetransmission rate of the upper layer will equal the transmission rate ofthe original stream, minus the combined rate of all lower layers.Because the transmission rate of the lower layers may have a fixedmaximum, the upper layer's rate may fluctuate and follow the variationsof the original stream.

In one embodiment, if the tokens all rotate at the same speed, this doesnot imply that they visit the same symbol slots at the same time.Instead, they will usually be at different positions in their ring. Thisis a result of empty slots that are skipped without waiting.

In one embodiment, the layered thinning algorithm has out of ordertransmission and may cause additional delay. Every ring can contain anupdate for the same symbol at a certain time and because the tokens canhave different column locations inside their ring, it is undefined whichring will dequeue and transmit its update first, which will be secondand so on. This means that not only the order of symbols is shuffled,but even the updates for a single symbol are no longer guaranteed to betransmitted chronologically.

While new updates for a symbol are queued, shifted and sent out, in somecases an update can survive in the rings substantially longer than itsyounger successors. This can occur when an update that is just about tobe dequeued by the token, gets shifted to the next layer by a morerecent update that entered through the lower ring. In this case it islikely that the more recent update is dequeued sooner than the olderupdate that now has to wait for the token of its new ring to reach itsslot. To address this issue, the system provides a way of putting theoutgoing, layered updates back in the original, chronological order.This is done by labeling incoming updates with sequential numbers andhaving the tokens pass their dequeued, labeled updates to a transmitwindow that uses the sequence numbers to restore the chronologicalorder. When updates are shifted out of the upper ring and thereforeconsidered lost, the transmit window is notified not to wait for theirsequence numbers.

When applying order reconstruction across all rings, the transmit windowmust have the same capacity as all rings combined, size_(nv)=N·j, wheresize_(nv) is the maximum capacity of the transmit window, N the numberof symbols in the original stream (and therefore the number of slots perring) and j the number of layers used. The potential size of thetransmit window, together with its bursty behavior when a delayed quoteupdate releases a substantial amount of pending updates from the window,may justify a mechanism that applies order reconstruction to achievechronologic transmission per symbol only, similar to the output of thebase layer encoder.

Like the algorithm depicted in FIG. 11A, the layered system introducesdelay. The model is essentially a collection of instances of the firstalgorithm, thus, its worst case delay is the sum of each ring's worstcase delay:

$t_{\max} = {\sum\limits_{i = 0}^{j}\frac{N}{f_{update}(i)}}$

where t_(max) is the maximum delay in seconds that the algorithm couldintroduce for a single quote update, N the number of symbols in theoriginal stream, j the number of used layers and f_(update)(i) theupdate frequency of the token in layer i.

In the foregoing specification, the invention has been described withreference to specific embodiments thereof. It will however, be evidentthat various modification and changes can be made thereto withoutdeparting from the broader spirit and scope of the invention as setforth in the appended claims. The specification and drawings are,accordingly, to be regarded in an illustrative rather than a restrictivesense.

What is claimed is:
 1. A video router, comprising: a memory; and aprocessor, wherein the processor executes instructions stored in thememory to cause the video router to: receive a layered video data streamincluding a base layer and a set of enhancement layers, identifybandwidth-limited conditions of an internet protocol network between thevideo router and a plurality of video receivers, forward the base layerfrom the video router to at least two of the plurality of videoreceivers via the internet protocol network, and selectively forward oneor more of the set of enhancement layers, but fewer than all of the setof enhancement layers, to at least two of the plurality of videoreceivers through the internet protocol network based upon theidentified bandwidth-limited conditions, and wherein the video routertransmits the layered video data stream according to an internetprotocol; wherein each layer of the layered video data stream comprisesdata packets, each of which is encoded with a sequence number and alayer identifier, and wherein the layer identifier for each data packetis based upon a layer to which the packet belongs.
 2. The video routerof claim 1, wherein the video router further selectively forwards theone or more of the set of enhancement layers based on video processingcapabilities of the video receiver.
 3. The video router of claim 1,wherein the video router divides available bandwidth among a pluralityof data streams, and wherein the video router prioritizes forwarding ofthe set of enhancement layers.
 4. The router of claim 1, wherein thevideo router further selectively forwards the one or more of the set ofenhancement layers based on a request for retransmission of one or morepackets from the layered video data stream.
 5. The video router of claim1, further comprising: a communications link connecting the video routerwith at least a second video router, wherein the layered video datastream is transmitted over the communications link between the videorouter and the second video router.
 6. A method for transmitting videosignals, the method comprising: receiving a layered video data streamcomprising a base layer and a set of enhancement layers; identifyingbandwidth-limited conditions of an internet protocol network between avideo router and a plurality of video receivers; forwarding the baselayer to at least two of the plurality of video receivers via theinternet protocol network; and selectively forwarding one or more of theset of enhancement layers, but fewer than all of the set of enhancementlayers, to at least two of the plurality of video receivers through theinternet protocol network based upon the identified bandwidth-limitedconditions; wherein the layered video data stream is transmittedaccording to an internet protocol; and wherein each layer of the layeredvideo data stream comprises data packets, each of which is encoded witha sequence number and a layer identifier, and wherein the layeridentifier for each data packet is based upon a layer to which thepacket belongs.
 7. The method of claim 6, wherein the selectivelyforwarding step further comprises: selectively forwarding the one ormore of the set of enhancement layers based on video processingcapabilities of the video receiver.
 8. The method of claim 6, furthercomprising: dividing available bandwidth among a plurality of datastreams; and prioritizing the set of enhancement layers for theselectively forwarding step.
 9. The method of claim 6, wherein theselectively forwarding step further comprises: selectively forwardingthe one or more of the set of enhancement layers based on a request forretransmission of one or more data packets from the layered video datastream.
 10. The method of claim 6, wherein the layered video data streamis received from another video router.
 11. A scalable video codingrouter, comprising: a memory; and a processor, wherein the processorexecutes instructions stored in the memory to cause the scalable videocoding router to: receive a layered video data stream that comprises abase layer and a set of enhancement layers, identify bandwidth-limitedconditions of an internet protocol network between the video router anda set of video receivers, and forward the base layer from the videorouter to at least two of the set of video receivers; wherein thescalable video coding router forwards all of the set of enhancementlayers to at least two of the video receivers in the set of videoreceivers with bandwidth-sufficient conditions; wherein the scalablevideo coding router selectively forwards one or more of the set ofenhancement layers, but fewer than all of the set of enhancement layers,to at least two of the remaining video receivers in the set of videoreceivers based upon the identified bandwidth-limited conditions;wherein the layered video data stream is transmitted according to aninternet protocol; and wherein each layer of the layered video datastream comprises data packets, each of which is encoded with a sequencenumber and a layer identifier, and wherein the layer identifier for eachdata packet is based upon a layer to which the packet belongs.
 12. Thevideo router of claim 11, wherein the video router further selectivelyforwards the one or more of the set of enhancement layers based on videoprocessing capabilities of each of the set of video receivers.
 13. Thevideo router of claim 11, wherein the video router divides availablebandwidth among a plurality of data streams, and wherein the videorouter prioritizes forwarding of the set of enhancement layers.
 14. Thevideo router of claim 11, wherein the video router further selectivelyforwards the one or more of the set of enhancement layers based on arequest for retransmission of one or more packets from the layered videodata stream.
 15. The video router of claim 11, further comprising: acommunications link connecting the video router with at least a secondvideo router, wherein the layered video data stream is transmitted overthe communications link between the video router and the second videorouter.